reperiendi

CCCs and lambda calculus

Posted in Category theory, Math, Programming, Quantum by Mike Stay on 2006 August 22

I finally got it through my thick skull what the connection is between lambda theories and cartesian closed categories.

The lambda theory of a set has a single type, with no extra terms or axioms. This gives rise to a CCC where the objects are generated by products and exponentials of a distinguished object, the free CCC on one object. For example, from a type X we generate objects 1, X, XX, (X^X)^{(X^X)} = X^{XX^X}, etc. The morphisms are reduction-equivalence classes of all the programs we can make out of the appropriately-typed S and K combinators. A product-exponential-preserving functor from this CCC to the category Set (which is also a CCC) picks out a set S_X for X, and because it preserves the structure, maps the product AB to S_A \times S_B and the exponential A^B to {S_A}^{S_B}.

The functor itself, however, can be uncomputable: one could, for example, have S_X be the set of halting programs for some universal Turing machine. This set is only computably enumerable, not computable.

When we have types and axioms involved, then we add structure to the set, constraints that the sets and functions on them have to satisfy. For instance, in the lambda theory of groups, we have:

  • a type G
  • terms
    • m:XX \to X
    • inv:X\to X
    • e:1 \to X
  • axioms for right- and left-unit laws, right-and left-inverses, and associativity

The CCC arising from this theory has all the morphisms from the free CCC on one object and extra morphisms arising from products and compositions of the terms. A structure-preserving functor to Set assigns G to a set and m, inv, and e to functions satisfying the axioms. These functions needn’t be computable, either—they only have to satisfy the group axioms.

So in terminology programmers are more familiar with, the terms and axioms define an abstract data type, an interface. The functor gives a class implementing the interface. But this implementation doesn’t need to be computable! Here’s another example: we start with a lambda theory with a data type \mathbb{N}, along with a term succ:\mathbb{N}\to \mathbb{N} and the axioms of Peano arithmetic; a functor from this lambda theory to Set will give us an implementation of natural numbers. Now we add a term f:\mathbb{N} \to \mathbb{N} to the theory, with no other constraints. One model of this theory is Peano arithmetic with an oracle to \Omega: it assigns f to the function that returns the nth bit of the Omega number for a universal prefix-free Turing machine.

I think that in order to get a computable model, we have to use a “computability functor” (my term). If I’m right, this means that instead of taking a functor directly into Set, we have to take a functor into a CCC with no extra terms to get a “computable theory” (again, my term), and then another from there into Set. This way, since all the morphisms in the category arising from the computable theory are built out of S and K combinators, the functor has to pick an explicit program for the implementation, not just an arbitrary function. From there, whether the implementation of the S and K combinators is computable or not really doesn’t matter; we can’t get at anything uncomputable from within the language.

Now, changing gears and looking at the “programs as proofs” aspect of all this: morphisms in the free CCC on one object are proofs in a minimal intuitionistic logic, where \to now means implication rather than exponentiation. The only two axioms are the ones from S and K. Adding a term of a given type to the theory adds a new axiom to the logic, while adding an axiom to the theory defines an equivalence of proofs in the logic.

A computable theory wouldn’t add any axioms, just assign names to proofs so they can be used as subproofs. And because the programs already satisfy the axioms of the computable theory, asserting the equivalence of two proofs is redundant: they’re already equivalent.

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